U.S. patent number 7,653,740 [Application Number 11/180,720] was granted by the patent office on 2010-01-26 for method and system for bandwidth allocation tracking in a packet data network.
This patent grant is currently assigned to Arris Group. Invention is credited to Adisak Mekkittikul, William J. Tuohy, Nader Vijeh.
United States Patent |
7,653,740 |
Mekkittikul , et
al. |
January 26, 2010 |
Method and system for bandwidth allocation tracking in a packet
data network
Abstract
In a metropolitan area network, a method and system for
maintaining an accurate total of the amount of allocated bandwidth
on the network. A plurality of incoming packets are assigned to a
respective plurality of queues of a metropolitan area network
switch. Using a fair arbitration scheme, the respective queues are
configured to empty at a specified output rate. A finish time for
each respective queue is computed, the finish time describing a
time at which the respective queue will be emptied using the output
rate. The plurality of queues are grouped into multiple groups in
accordance with their respective finish times. The earliest group
includes the reserved rates of those queues having a finish time
indicating an empty condition at a first time increment. The second
earliest group includes the reserved rates of those queues having a
finish time indicating an empty condition at a second time
increment later than the first time increment, and so on. The
amount of allocated bandwidth on the network is determined by
tracking the sum of the reserved rates of all the multiple groups.
The first time increment, second time increment, and the like are
indexed with respect to a schedule clock. The earliest group thus
indicates those queues that will have an empty condition at a next
time increment of the schedule clock. The determination of the
amount of allocated bandwidth can be accomplished in real time,
thereby allowing the efficient allocation of unallocated bandwidth
in real time.
Inventors: |
Mekkittikul; Adisak (Mountain
View, CA), Vijeh; Nader (Sunnyvale, CA), Tuohy; William
J. (Sunnyvale, CA) |
Assignee: |
Arris Group (Suwanee,
GA)
|
Family
ID: |
34067621 |
Appl.
No.: |
11/180,720 |
Filed: |
July 14, 2005 |
Prior Publication Data
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Document
Identifier |
Publication Date |
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US 20050249128 A1 |
Nov 10, 2005 |
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Related U.S. Patent Documents
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Application
Number |
Filing Date |
Patent Number |
Issue Date |
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10094035 |
Mar 7, 2002 |
6947998 |
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60274621 |
Mar 8, 2001 |
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Current U.S.
Class: |
709/235; 709/240;
370/412 |
Current CPC
Class: |
H04L
12/2852 (20130101); H04L 47/2441 (20130101); H04L
47/6255 (20130101); H04L 12/42 (20130101); H04L
47/522 (20130101); H04L 47/30 (20130101); H04L
47/50 (20130101); H04L 47/562 (20130101); H04L
47/2433 (20130101); H04L 47/6215 (20130101) |
Current International
Class: |
G06F
15/16 (20060101); H04L 12/28 (20060101) |
Field of
Search: |
;709/214,240,242,235
;370/468,477,232-235,253,412 |
References Cited
[Referenced By]
U.S. Patent Documents
Primary Examiner: Lim; Krisna
Attorney, Agent or Firm: FSP LLC
Parent Case Text
This application claims the benefit of earlier filed U.S.
Provisional Application "A METHOD AND SYSTEM FOR BANDWIDTH
ALLOCATION TRACKING IN AN ASYNCHRONOUS METRO PACKET TRANSPORT RING
NETWORK", by Mekkittikul et al., Ser. No. 60/274,621, filed on Mar.
8, 2001
Claims
What is claimed is:
1. A network device comprising: logic to group individual data
flows into buckets, each bucket comprising a plurality of queues;
and logic to track an active and inactive status of the individual
data flows in real time.
2. The network device of claim 1, wherein the logic to group
individual data flows into buckets, each bucket comprising a
plurality of queues further comprises: each bucket comprising
queues having a same empty time.
3. The network device of claim 1, further comprising: logic to
allocate the bandwidth of inactive flows to active flows in real
time.
4. The network device of claim 3, wherein the logic to allocate the
bandwidth of inactive flows to active flows in real time further
comprises: logic to allocate the bandwidth based at least in part
on finish times of the queues.
5. The network device of claim 1, further comprising: logic to
track a total allocated bandwidth for all active data flows.
6. A network comprising: one or more devices together comprising
logic to group individual data flows into buckets, each bucket
comprising a plurality of queues; and logic to track an active and
inactive status of the individual data flows in real time.
7. The network of claim 6, wherein the logic to group individual
data flows into buckets, each bucket comprising a plurality of
queues further comprises: each bucket comprising queues having a
same empty time.
8. The network of claim 6, further comprising: logic to allocate
the bandwidth of inactive flows to active flows in real time.
9. The network of claim 8, wherein the logic to allocate the
bandwidth of inactive flows to active flows in real time further
comprises: logic to allocate the bandwidth based at least in part
on finish times of the queues.
10. The network of claim 6, further comprising: logic to track a
total allocated bandwidth for all active data flows.
11. A process in a network device, comprising: grouping individual
network data flows into buckets implemented as data storage
locations within the network device, each bucket comprising a
plurality of queues; and the network device tracking an active and
inactive status of the individual data flows in real time.
12. The process of claim 11, wherein grouping individual data flows
into buckets, each bucket comprising a plurality of queues further
comprises: grouping the flows into buckets each bucket comprising
queues having a same empty time.
13. The process of claim 11, further comprising: allocating the
bandwidth of inactive flows to active flows in real time.
14. The process of claim 13, wherein allocating the bandwidth of
inactive flows to active flows in real time further comprises:
allocating the bandwidth based at least in part on finish times of
the queues.
15. The process of claim 11, further comprising: tracking a total
allocated bandwidth for all active data flows.
Description
FIELD OF THE INVENTION
The present invention relates to the field of packet data networks.
More specifically, the present invention pertains a data flow
control method and system for managing the data flow with respect
to the available bandwidth in a metro packet transport ring
network.
BACKGROUND ART
The Internet is a general purpose, public computer network which
allows millions of computers all over the world, connected to the
Internet, to communicate and exchange digital data with other
computers also coupled to the Internet. As new technologies emerge,
the speed at which one can connect onto the Internet is ever
increasing. Now, users on the Internet have the bandwidth to
participate in live discussions in chat rooms, play games in
real-time, watch streaming video, listen to music, shop and trade
on-line, etc. In the future, it is imagined that the bandwidth will
be such that video-on-demand, HDTV, IP telephony, video
teleconferencing, and other types of bandwidth intensive
applications will soon be possible.
One approach by which bandwidth is being increased relates to fiber
optics technology. By sending pulses of light through glass fibers
no thicker than a human hair, vast amounts of digital data can be
transmitted at extremely high speeds. And with the advent of dense
wavelength division multiplexing, different wavelengths of light
can be channeled over the same, single fiber strand, thereby
increasing its capacity several fold.
However, there is a problem with distributing the bandwidth of this
new fiber optic network to end users. As is well known, some
applications are insensitive to bandwidth constraints, such as
latency and dropped packets. For example, email applications and
basic Web browsing are relatively time insensitive to latency and
dropped packets. On the other hand, applications such as real time
two way voice communication or video are very sensitive to time
delays caused by latency and dropped packets. Acceptable
performance of these applications is highly dependent upon the
provision of a guaranteed minimum bandwidth.
Unfortunately, due to network traffic congestion, network
availability, routing conditions, and other uncontrollable external
factors, the provisioning of a guaranteed level of bandwidth
availability for certain customers has proven problematic. In
general, data packets vie for available bandwidth and are routed
according to a best-effort delivery model. As such, the reliability
of traditional packet switched data networks is at times
sub-optimal. For example, in most cases, it is very difficult to
provide any kind of quality of service (QoS) using traditional LAN
switches and routers on IP networks. QoS refers to the guarantee of
providing timely delivery of information, controlling bandwidth per
user, and setting priorities for select traffic.
Different network traffic flows (or simply "flows") are
respectively associated with different applications. A flow refers
to the transmission of packets from a sender to a receiver to
support an application, such as transferring a Web page,
implementing a voice over IP conversation, playing a video, or the
like. Some flows are described as real time flows since they
require very low latency (e.g., a voice over IP application). Other
flows are not so much latency dependent as they are consistent data
transfer rate dependent (e.g., video over the Web). For real-time
application flows such as video on demand, HDTV, voice
communications, etc., dropped packets or late-arriving packets of
the flows can seriously disrupt or even destroy performance. And
for many Internet Service Providers (ISP's), Applications Service
Providers (ASP's), web sites/portals, and businesses, it is of
paramount importance that they have the ability to provide these
flows with a certain minimum threshold bandwidth and/or latency.
For example, an e-commerce or business web site may lose critical
revenue from lost sales due to customers not being able to access
their site during peak hours.
Because QoS is so highly desired by some users, there are
mechanisms which have been developed to provide QoS functionality.
One prior art method for implementing QoS is the use of various TDM
(time division multiplexing) schemes. One widely used TDM scheme is
the implementation of T-carrier services (e.g., T1 line for
carrying data at 1.544 Mbits/sec. and T3 line for carrying data at
a much faster rate of 274.176 Mbits/sec). These T1 and T3 lines are
dedicated point-to-point datalinks leased out by the telephone
companies. The telephone companies typically charge long distance
rates (e.g., $1,500-$20,000 per month) for leasing out a plain old
T1 line. Another commonly used TDM scheme for achieving QoS relates
to Synchronous Optical Network (SONET). As with T-carrier services,
SONET uses TDM to assign individual channels, or flows, to
pre-determined time slots. With TDM, each channel is guaranteed its
own specific time slot in which it can transmit its data. Although
TDM enables QoS, it is costly to implement because both the
transmitter and receiver must be synchronized at all times. The
circuits and overhead associated with maintaining this precise
synchronization is costly. Furthermore, TDM based networking
technologies are inefficient with respect to unused time slots. If
flows are inactive, their allocated bandwidth is wasted. In
general, with TDM technologies, unused bandwidth from inactive
flows is not reallocated to other users.
Another prior art method is the use of various forms of bandwidth
reservations in conjunction with asynchronous schemes. Asynchronous
data transmission schemes provide numerous advantages when compared
to synchronous TDM type schemes, and as such, are generally
overtaking synchronous technologies in both voice and data network
installations (e.g., the IP based networks of the Internet). In
implementing QoS, asynchronous schemes usually function by
reserving a portion of their bandwidth for "high priority" latency
sensitive flows. With most asynchronous schemes (e.g., Ethernet),
QoS performance deteriorates with the increasing bandwidth
utilization of the network. As the percentage of available
bandwidth utilized by the network increases, the less efficient the
prior art asynchronous QoS reservation schemes perform. Such
schemes either maintain a large margin of unused bandwidth to
ensure QoS, thereby virtually guaranteeing an under-utilization of
available total bandwidth, or over-allocate bandwidth, leading to
the abrupt dropping data for some users and/or ruining QoS for any
high priority users.
Thus what is required is a solution that provides the advantages of
asynchronous data networks while efficiently implementing QoS. What
is required is a solution that enables the efficient allocation of
available bandwidth, thereby allowing guaranteed QoS. The required
solution should be able to allocate bandwidth to individual flows
asynchronously without incurring the deteriorating performance of
prior art asynchronous schemes with increasing scale (e.g.,
extremely large number of flows) and increasing network
utilization. The required solution should be able to ensure a
minimum amount of reserved bandwidth without incurring the wasted
bandwidth problems of prior art TDM based networking schemes where
bandwidth is wasted on inactive flows.
The required solution should be able to track individual flows on
an individual basis, in order to ensure individual flows are not
starved of bandwidth, while simultaneously ensuring bandwidth is
not over-allocated to flows which do not require it. The required
solution should be able to track when individual flows are active
and when they are inactive, thereby allowing the bandwidth
allocated to the inactive flows to be reassigned to those flows in
need of it. The required solution should be capable of tracking
total allocated bandwidth in real time, thereby allowing efficient
allocation of unused bandwidth in real time while maintaining QoS.
The real-time total allocated bandwidth tracking should allow the
dynamic allocation of unused bandwidth in real-time. The present
invention provides a novel solution to the above requirements.
SUMMARY OF THE INVENTION
The present invention comprises a method and system that provides
the advantages of asynchronous data networks while efficiently
implementing QoS. The present invention enables the efficient
allocation of available bandwidth, thereby allowing guaranteed QoS.
The present invention is able to track individual flows on an
individual basis, in order to ensure individual flows are not
starved of bandwidth, while simultaneously ensuring bandwidth is
not over-allocated to flows which do not require it. The present
invention can track when individual flows are active and when they
are inactive, thereby allowing the bandwidth allocated to the
inactive flows to be reassigned to those flows in need of it. The
present invention can track total allocated bandwidth in real time,
thereby allowing efficient allocation of unused bandwidth in real
time while maintaining QoS. The real-time total allocated bandwidth
tracking allows the dynamic allocation of unused bandwidth in
real-time, while maintaining QoS.
In one embodiment, the present invention is a system for
maintaining an accurate total of the amount of allocated bandwidth
on the network, as implemented within a metropolitan area switch
(MPS) that functions by allocating bandwidth of a metropolitan area
network. Within the MPS, a plurality of incoming packets are
assigned to a respective plurality of queues of the MPS. A finish
time for each respective queue is computed, the finish time
describing a time at which the respective queue will be emptied
using the output rate. The plurality of queues are grouped into
multiple groups in accordance with their respective finish times.
These groups are referred to as "buckets" due to the fact that they
include those queues having the same finish times.
The earliest group includes the reserved bandwidth of those queues
having a finish time indicating an empty condition at a first time
increment. The second earliest group includes the reserved
bandwidth of those queues having a finish time indicating an empty
condition at a second time increment later than the first time
increment, and so on. Thus, for example, bucket 0 contains those
queues which will be empty at the next time increment, bucket 1
contains those queues that will be empty at the next two time
increments, and so on. The amount of allocated bandwidth on the
network is determined by counting the reserved bandwidth of all
active flows.
The first time increment, second time increment, and the like are
indexed with respect to a schedule clock. One increment of the
schedule clock comprises one complete round robin arbitration
(e.g., per queue output onto the metropolitan area network) of all
active queues within the MPS. The earliest group thus indicates
those queues that will have an empty condition at a next time
increment (e.g., output round) of the schedule clock. A new finish
time is computed for each respective queue when a new packet is
received by the respective queue. In this manner, the series of
buckets are progressively "emptied" as the schedule clock
progresses, and new buckets are filled as new queues receive new
packets for transmission and new associated empty times. The queues
that are empty at the next time increment indicate those flows that
will be inactive at the next time increment. The bandwidth
allocated to those flows can be reallocated. In this manner, the
determination of the amount of allocated bandwidth can be
accomplished in real time, thereby allowing the efficient
allocation of unallocated bandwidth in real time while maintaining
quality of service. The earliest bucket (e.g., bucket 0) shows the
reserved rate of all queues which will be empty in the next time
increment.
Thus, by grouping individual flows into buckets as described above,
embodiments of the present invention can efficiently scale up to
handle an extremely large number (e.g., 1 million or more)
individual flows. The flows are assigned to buckets as described
above on an individual basis. Their condition (active vs. inactive)
is individually tracked in real-time, allowing their allocated
bandwidth for inactive flows to be reallocated to active flows in
real time. In so doing, the present invention enables the efficient
allocation of available bandwidth, since the MPS is capable of
tracking total allocated bandwidth in real time. This allows the
efficient allocation of unused bandwidth in real time while
maintaining QoS.
BRIEF DESCRIPTION OF THE DRAWINGS
The present invention is illustrated by way of example and not by
way of limitation, in the figures of the accompanying drawings and
in which like reference numerals refer to similar elements and in
which:
FIG. 1 shows the overall architecture of the asynchronous metro
packet transport ring network according to the currently preferred
embodiment of the present invention.
FIG. 2 shows an exemplary Metro Packet Transport Ring.
FIG. 3 shows an exemplary diagram of components of an MPTR.
FIG. 4 a diagram of a set of MPS units and ring segments as
implemented within an exemplary system in accordance with one
embodiment of the present invention.
FIG. 5 shows a diagram of a queue of an MPS and its associated
finish time.
FIG. 6A shows a diagram depicting the multi-group queuing process
in accordance with one embodiment of the present invention.
FIG. 6B graphically depicts the summation of all r.sub.i and
w.sub.i in accordance with one embodiment of the present
invention.
FIG. 7 shows a diagram of a bucket information base (BIB) in
accordance with one embodiment of the present invention.
FIG. 8 shows a flow information base (FIB) in accordance with one
embodiment of the present invention.
FIG. 9 shows a flow chart of the steps of a bandwidth tracking and
allocation process in accordance with one embodiment of the present
invention.
DETAILED DESCRIPTION OF THE INVENTION
Reference will now be made in detail to the embodiments of the
invention, examples of which are illustrated in the accompanying
drawings. While the invention will be described in conjunction with
the preferred embodiments, it will be understood that they are not
intended to limit the invention to these embodiments. On the
contrary, the invention is intended to cover alternatives,
modifications and equivalents, which may be included within the
spirit and scope of the invention as defined by the appended
claims. Furthermore, in the following detailed description of the
present invention, numerous specific details are set forth in order
to provide a thorough understanding of the present invention.
However, it will be obvious to one of ordinary skill in the art
that the present invention may be practiced without these specific
details. In other instances, well known methods, procedures,
components, and circuits have not been described in detail as not
to unnecessarily obscure aspects of the present invention.
Embodiments of the present invention are directed to a method and
system for maintaining an accurate total of the amount of allocated
bandwidth on a network, as implemented within a metropolitan area
switch (MPS). The present invention provides the advantages of
asynchronous data networks while efficiently implementing QoS. The
present invention enables the efficient allocation of available
bandwidth, thereby allowing guaranteed QoS. The present invention
is capable of tracking total allocated bandwidth in real time,
thereby allowing efficient allocation of unused bandwidth in real
time while maintaining QoS. The present invention and its benefits
are further described below.
FIG. 1 shows an overall architecture of an asynchronous metro
packet transport ring network in accordance with a currently
preferred embodiment of the present invention. A metropolitan
packet transport ring (MPTR) consists of a ring which is laid to
transmit data packets in a metropolitan area network (MAN). A MAN
is a backbone network which spans a geographical metropolitan area.
Typically, telephone companies, cable companies, and other
telecommunications providers supply MAN services to other
companies, businesses, and users who need access to networks
spanning public rights-of-way in metropolitan areas. In the
currently preferred embodiment, the communications channel of the
MPTR is implemented using a ring topology of installed fiber optic
cables. Other less efficient transmission mediums such as hybrid
fiber coax, coax cables, copper wiring, or even wireless (radio
frequency or over-the-air laser beams) can be used or substituted
in part thereof.
Users coupled to a particular MPTR can transmit and receive
packetized data to/from each other through that MPTR. For example,
a personal computer coupled to MPTR1 can transmit and received data
packets to/from a server also coupled to MPTR1. Furthermore, data
packets originating from one MPTR can be routed to another MPTR by
means of a router. For example, a computer coupled to MPTR1 can
transmit data packets over its fiber ring to a router 101 to MPTR2.
The data packets can then be sent to its final destination (e.g., a
computer coupled to MPTR2) through the fiber ring associated with
MPTR2. It should be noted that the MPTR rings can be of various
sizes and configurations. Although the currently preferred
embodiment contemplates the use of a ring, the present invention
can also utilize other types of topologies. The MPTRs can also be
coupled onto the Internet backbone via a router. For example, MPTR1
can be coupled to a dense wavelength division multiplexed (DWDM)
fiber backbone 102 by means of router 101. Thereby, users coupled
to MPTR1 has access to the resources available on traditional
Internet 103. Note that the present invention can be used in
conjunction with traditional Internet schemes employing standard
routers, switches, and other LAN equipment 104-107. And any number
of MPTR's can thusly be coupled together to gracefully and
cost-efficiently scale to meet the most stringent networking
demands which may arise. And as one particular ring becomes
overloaded, a second, third, forth, etc. MPTR may be added to
accommodate the increased load. These MPTR's can be coupled to the
same router (e.g., MPTR5, MPTR6, and MPTR7) or may alternatively be
coupled to different routers.
Not only does the present architecture scale gracefully, it also
offers great flexibility. In one embodiment, an MPTR can be used to
support one or more LANs. For instance, MPTR6 may support traffic
flowing to/from LAN 108. Optionally, it is conceivable that an MPTR
may be coupled directly to another MPTR. In this manner, data
flowing in MPTR8 can be directly exchanged with data packets
flowing through MPTR7. Alternatively, a single MPTR can have
multiple entries/exits. For example, MPTR5 is coupled to both
router 109 as well as router/switch 110. Thereby, users on MPTR5
have the ability to transmit and receive data packets through
either of the two routers 109 or 110. Virtually any configuration,
protocol, medium, and topology is made possible with the present
MPTR invention.
The implementation and functionality of an MPTR is now described.
Referring to FIG. 2, an exemplary Metro Packet Transport Ring 200
is shown. It can be seen that MPTR 200 is comprised of two fiber
cable rings, or rings, 201 and 202; a number of Metro Packet
Switches (MPS1-MPSn); and a Ring Management System (RMS) 203. The
physical layer of an MPTR is actually comprised of two redundant
fiber cable rings 201 and 202. Data packets flow in opposite
directions through the two rings (e.g., clockwise in ring 201 and
counter-clockwise in ring 202). Dispersed along the fiber rings 201
and 202 are a number of Metro Packet Switches (MPS's). An MPS is
coupled to both of the fiber rings 201 and 202. Thereby, if there
is a break in one segment of the fiber ring, data can be redirected
through one of the MPS's to flow through the other, operational
fiber ring. Alternatively, traffic can be re-directed to minimize
localized congestion occurring in either of the rings.
In the currently preferred embodiment, each MPTR can support up to
254 MPS's. An MPS is a piece of equipment which can be housed in
specially designed environmental structures or it can be located in
wiring closets or it can reside at a place of business, etc. The
distances between MPS's can be variable. It is through an MPS that
each individual end user gains access to the fiber rings 201 and
202. Each individual end user transmits packetized data onto the
MPS first. The MPS then schedules how that packetized data is put
on the fiber ring. Likewise, packetized data are first pulled off a
fiber ring by the MPS before being sent to the recipient end user
coupled to the MPS. In the currently preferred embodiment, a single
MPS can support up to 128 end users. An end user can be added to an
MPS by inserting a line interface card into that particular MPS.
The line interface cards provide I/O ports through which data can
be transferred between the MPS and its end users. Different line
interface cards are designed in order to meet the particular
protocol corresponding to that particular end user. Some of the
protocols supported include T1, T3, SONET, Asynchronous Transfer
Mode (ATM), digital subscriber line (DSL) Ethernet, etc. It should
be noted that line interface cards can be designed to meet the
specifications of future protocols. In this manner, end users such
as mainframe computers, workstations, servers, personal computers,
set-top boxes, terminals, digital appliances, TV consoles, routers,
switches, hubs, and other computing/processing devices, can gain
access to either of the fiber rings 201 and 202 through an MPS.
Not only does an MPS provide I/O ports to end users, but an MPS
also provides a means for inputting packetized data into the MPTR
and also for outputting packetized data out from the MPTR. For
example, data packets are input to MPTR 200 via MPS 204 which is
coupled to router 205. Similarly, data packets are output from MPTR
200 via MPS 204 to router 205.
Another function of an MPS entails passing along incoming data
packets originating from an upstream MPS to the next downstream
MPS. An MPS receives upstream data packets forwarded from an
upstream MPS via an input fiber port coupled to the fiber ring.
Data packets received from the fiber ring are examined by that MPS.
If the data packet is destined for an end user coupled to that
particular MPS, the data packet is routed to the appropriate I/O
port. Otherwise, the MPS immediately forwards that data packet to
the next downstream MPS as quickly as possible. The data packet is
output from the MPS by an output fiber port onto the fiber ring. It
should be noted that such pass-through packets flowing from an
upstream fiber ring segment, through the MPS, and onto a downstream
fiber ring segment, always takes priority over packets waiting to
be inserted onto the fiber ring by the MPS. In other words, the MPS
inserts data packets generated by its end users only as bandwidth
permits.
An example is now offered to show how data packets flow in an MPTR.
With reference to FIG. 2, a computer 207 coupled to MPS 4 can
transmit and receive data to/from the Internet as follows. Data
packets generated by the computer are first transmitted to MPS 4
via a line coupled to a line interface card residing within MPS4.
These data packets are then sent on to MPS3 by MPS4 via ring
segment 206. MPS3 examines the data packets and passes the data
packets downstream to MPS2 via ring segment 207; MPS2 examines the
data packets and passes the data packets downstream to MPS1 via
ring segment 208. Based on the addresses contained in the data
packets, MPS1 knows to output theses data packets on to the I/O
port corresponding to router 205. It can be seen that MPS1 is
connected to a router 205. Router 205 routes data packets to/from
MPTR 200, other MPTR's, and the Internet backbone. In this case,
the data packets are then routed over the Internet to their final
destination. Similarly, data packets from the Internet are routed
by router 205 to MPTR 200 via MPS1. The incoming data packets are
then examined and forwarded from MPS1 to MPS2 via ring segment 209;
examined and forwarded from MPS2 to MPS3 via ring segment 210; and
examined and forwarded from MPS3 to MPS4 via ring segment 211. MPS4
examines these data packets and determines that they are destined
for computer 207, whereby MPS4 outputs the data packets through its
I/O port corresponding to computer 207.
Likewise, users coupled to any of the MPS's can transmit and
receive packets from any other MPS on the same MPTR without having
to leave the ring. For instance, a user on MPS2 can transmit data
packets to a user on MPS4 by first transmitting the packets into
MPS2; sending the packets from MPS2 to MPS3 over ring segment 207;
MPS3 sending the packets to MPS4 over ring 202; and MPS4 outputting
them on the appropriate port corresponding to the intended
recipient.
Referring still to FIG. 2, it should be noted that the present
invention solves the strict priority problems common to ring
topology networks. The strict priority problem refers to the fact
that upstream nodes (e.g., an upstream MPS) have larger amounts of
available bandwidth in the communications channel in comparison to
downstream nodes. For example, in the case of ring segment 210, MPS
2 is able to insert its local input flows (e.g., insertion traffic)
onto segment 210 prior to MPS 3, and so on with MPS 3 and MPS 4
with ring segment 211. Hence, MPS 4, by virtue of its location
within the ring topology, has less available bandwidth to insert
its local input flow in comparison to MPS 3 and MPS 2.
To avoid strict priority problems, detailed information regarding
the allocated bandwidth of the ring segments is required. Each MPS
needs to be aware of the allocated bandwidth of the segments in
order to make intelligent decisions regarding the allocation of any
remaining unallocated bandwidth. Such information is even more
important where bandwidth utilization is to be maximized in
conjunction with guaranteed QoS. Preferably, bandwidth utilization
information should be available on a "per-flow" basis should be
sufficiently timely to allow intelligent allocation decisions to be
made in real time.
Additional descriptions of the architecture of the MPTR, MPS, and
RMS can be found in U.S. patent applications "GUARANTEED QUALITY OF
SERVICE IN AN ASYNCHRONOUS METRO PACKET TRANSPORT RING", filed on
Jun. 30, 2000, Ser. No. 09/608,747, assigned to the assignee of the
present invention which is incorporated herein in its entirety, and
"PER-FLOW CONTROL FOR AN ASYNCHRONOUS METRO PACKET TRANSPORT RING",
filed on Jun. 30, 2000, Ser. No. 09/608,489, assigned to the
assignee of the present invention which is incorporated herein in
its entirety.
FIG. 3 shows an exemplary diagram of components of an MPTR. A
number of MPS's 301-306 are shown coupled to a fiber ring 307. Two
of the MPS's 302 and 303 have been shown in greater detail to
depict how data flows in an MPTR. A number of computers 308-310 are
shown coupled to MPS 302. Each of these computers 308-310 has a
corresponding buffer 311-313. These buffers 311-313 are used to
temporarily store incoming data packets from their respective
computers 308-310. Associated with each of these buffers 311-313 is
a respective controller 314-316 which controls when packets queued
in that particular buffer are allowed to be transmitted onto the
ring 307. Once a packet is allowed to be transmitted out from MPS
302, it is inserted into an inserter 325 and added with the other
outbound packets for that cycle. Once a packet is conveyed from an
MPS onto ring 307, that packet is transmitted to its destination at
the maximum rate of ring 307 and immediately forwarded through
intermediary MPS's (if any).
In a preferred MPTR embodiment, fair bandwidth allocation is
implemented using a per-flow bandwidth allocation concept. Traffic
on the ring 307 is classified into flows. For example, all packets
from one user belong to one flow. The granularity of flow can be
fine (e.g., per-session) or coarse (e.g., per service port, etc.),
and is generally specifiable by packet classification rules. Once
packets are classified into a flow, each MPS can allocation
bandwidth to each flow fairly and monitor that no flow exceeds the
allocation.
The flow thus must be set up before the packets can be sent on the
ring. Setting up flow involves specifying a number of parameters.
Among these, the reserved bandwidth, r.sub.i, and the allocation
weight, w.sub.i, are necessary for flow control, where "i" is the
flow's unique identifier referred to as the flow ID. Once set up, a
flow is recognized by its unique flow ID.
FIG. 4 shows a diagram showing three MPS units and their respective
ring segments. As depicted in FIG. 4, three MPS units (MPS 0, MPS
1, and MPS 2) are shown with their respective ring segments
401-404. The MPS units are shown with their respective insertion
traffic (I0, I1, and I2) and their respective exit traffic (E0, E1,
and E2). Each MPS 0-2 is shown with a plurality of internal queues
(four depicted within each MPS) used for tracking the flows.
As shown in FIG. 4, the queues of each MPS tracks the allocated
bandwidth on each outgoing ring segment 401-404. As shown in FIG.
4, the traffic on the outgoing segment is represented as:
.times..times..times..times..times..times. ##EQU00001##
The queues of each MPS track the data traffic belonging to each
individual flow (described in greater detail below). The traffic on
each segment takes into account the exit traffic of the previous
MPS, the insertion traffic of the previous MPS, and the through
traffic on the ring. The insertion traffic of each MPS is shown in
FIG. 4 as "I" and the exit traffic of each MPS is shown as "E". The
insertion traffic is the flows from the users coupled to the MPS
that want to get onto the ring, for example, destined for users
coupled to another MPS. The exit traffic is the flows destined for
the users coupled to the MPS coming from other MPS units. The
queues within each MPS are used to track the unique flows (e.g.,
having unique flow IDs) that are monitored and maintained by an
MPS. Each of the queues tracking the outgoing flow for the outgoing
ring segment are drained at a rate equal to the allocated
bandwidth.
The queues are emptied at a rate affected by their respective
weight, w.sub.i. The weight of each queue allows the implementation
of differing levels of bandwidth per queue. For example, where
queues are of equal weight, the individual flow packets are routed
from the queues at an equal rate. Once a packet is inserted onto an
outbound ring segment, such as, for example, a packet from a flow
of insertion traffic I.sub.0 being inserted onto ring segment 402,
that packet is added with other outbound packets and is transmitted
along ring segment 402 at wire speed, or the maximum rate of
transmission of the ring. The packet is immediately forwarded
through intermediary MPS's (if any) as through traffic. Once a
queue becomes empty, its bandwidth allocation becomes available for
reassignment to other non-empty queues.
It should be noted that in a preferred embodiment, an MPS in
accordance with the present invention maintains large sets of
virtual queues (VQs) to monitor flow activity on all of its output
links. Virtual queues function in a manner similar to the queues
described above (e.g., the queues shown within the MPS units
depicted in FIG. 4), however, they are implemented as counters
which track the depth of the queues so that the data packets are
not delayed as flow through their respective buffers. Additional
descriptions of virtual queues as implemented in the preferred
embodiment can be found in "PER-FLOW CONTROL FOR AN ASYNCHRONOUS
METRO PACKET TRANSPORT RING", filed on May 13, 2004, publication
number US20050002392, which is incorporated herein in its entirety.
A VQ will have a finish time describing the time when all the
packets are completely drained from the VQ at a flow allocated rate
of f(subscript)i .
FIG. 5 shows a diagram of a queue 415 and its associated finish
time. The output rate of the queues 411-415 allows the
determination of a "finish time" describing the time at which the
respective queue will be emptied. This finish time provides a key
measure of the total allocated bandwidth of the ring 450. Thus, as
depicted in FIG. 5, queue 415 has a finish time that describes the
time at which queue 415 will be emptied at its output rate. When a
new packet arrives as shown, a new finish time is computed
reflecting the new depth of the queue 415. Thus, as depicted in
FIGS. 5 and 6, the MPS routes packets from the respective queues at
a specified output rate, and a finish time for each respective
queue is computed, the finish time describing a time at which the
respective queue will be emptied using the allocated output rate
(e.g., f.sub.i as defined below).
In this manner, each MPS maintains a large number of queues (e.g.,
up to 1 million or more), one for each flow at each link. Each
queue grows at the rate of the traffic belonging to the flow, and
is drained at a rate equal to the allocated bandwidth. Congestion
is measured in the forms of: .SIGMA.r.sub.i and .SIGMA.w.sub.i of
all non-empty (active) queues (e.g., queues 411-415). High values
of
.times..times..times..times..times..times. ##EQU00002## indicate
that more flows are competing for the outgoing link bandwidth of
the MPS. Each MPS frequently monitors the states of its queues to
update these two parameters. Once detected, an MPS uses
.times..times..times..times..times..times. ##EQU00003## to
calculate bandwidth allocation for each flow.
In a preferred embodiment, each MPS calculates a fair allocation of
bandwidth for all flows going through each of congestion points
(e.g., at the outgoing ring segments). The allocation is calculated
based on the following calculation:
.function..times. ##EQU00004## where f.sub.i denotes the allocated
bandwidth for flow i, and C is the link capacity of the congested
point. Note that the term
.times. ##EQU00005## is simply the unreserved bandwidth portion of
the link that the MPS needs to reallocate fairly based on the
reserved weights. This term is graphically depicted in FIG. 6B
below.
For bandwidth efficiency, each MPS does not send out f.sub.i for
every flow it sees. Instead, it sends a capacity reserved ratio
(CRR) which generally describes the amount of unallocated bandwidth
of the link. The CRR can then be used by each source within each
MPS to calculate its own f.sub.i from its static database of
r.sub.i and w.sub.i. CRR is more formally defined as follows:
.times. ##EQU00006## CRR is broadcasted to all other MPSs
periodically to enable all MPSs to allocate unallocated link
bandwidth. Each MPS can independently choose the frequency of the
update. For each received CRR, each source uses the equation below
to calculate its f.sub.i. f.sub.i=r.sub.i+w.sub.i*CRR
Thus, in order to efficiently distribute unallocated link
bandwidth, each MPS needs to track the total amount of allocated
bandwidth and the total weight of the allocated bandwidth,
.times..times..times..times..times..times. ##EQU00007## In
accordance with the present invention, these terms are tracked in
real time and track flow activity at high speeds, as high as 10
Gbps per ring segment. The present invention uses the finish times
of the respective queues and the assigned weights of the respective
queues to implement a high speed tracking method for
.times..times..times..times..times..times. ##EQU00008## These
techniques involve the uses of per-flow queues, a flow information
base (FIB), a bucket information base (BIB), and a schedule clock.
Using these terms, embodiments of the present invention can
efficiently scale up to handle an extremely large number (e.g., 1
million or more) individual flows, while remaining within the
capabilities of integrated circuit technology (e.g., can be
implemented in an ASIC). The individual flows can be tracked in
real-time, allowing their allocated bandwidth for inactive flows to
be reallocated to active flows in real time.
Referring now to FIG. 6A, a diagram depicting the multi-group
queuing process of the present embodiment is shown. FIG. 6A depicts
a plurality of flows sorted into a plurality of groups, shown as
bucket 0, bucket 1, bucket 2, and so on, to bucket n. The plurality
of queues are grouped into the multiple buckets, or groups, in
accordance with their respective finish times. The finish times are
indexed with respect to a schedule clock. The schedule clock, or
global clock, provides the time reference for finish times. The
value of schedule clock represents the current virtual time that
finish times are compared to. Schedule clocks increment at a rate
proportional to the congestion at a node, as described below. As
depicted in FIG. 6A, as buckets are emptied, they move from right
to left, as each bucket successively reaches the "queue empty"
state shown on the left side FIG. 6A.
These groups of flows are referred to as "buckets" due to the fact
that they include those queues having the same finish times with
the schedule clock. For example, bucket 0 includes the reserved
bandwidth and weight of those flows having a finish time
corresponding to the next increment of the schedule clock, while
bucket n includes the reserved bandwidth and weight of those flows
having the longest finish time with respect to the schedule clock.
Thus, the earliest bucket (e.g., bucket 0) includes those flows
(e.g., queues) having a finish time indicating an empty condition
at a first time increment, the second earliest bucket (e.g., bucket
1) includes those queues having a finish time indicating an empty
condition at a second time increment later than the first time
increment, and so on. Thus, for example, bucket 0 contains the
reserved bandwidth and weight of those queues which will be empty
at the next time increment of the schedule clock, bucket 1 contains
those queues that will be empty at the next two time increments of
the schedule clock, and so on, thereby indicating the amount of
unallocated bandwidth that becomes available each time increment.
The amount of allocated bandwidth on the network is determined by
counting the total allocated bandwidth and total allocated weight
of all the active flows (e.g., all bucket totals).
The time increments for the first bucket, the second bucket, and
the like are indexed with respect to the schedule clock. One
increment of the schedule clock comprises one complete round robin
arbitration (e.g., per queue output onto the metropolitan area
network) of all active queues within the MPS, in the manner
described in FIG. 4 above. Inactive, or empty, queues do not
contribute to the schedule clock period. The bucket 0 thus
indicates those flows that will have an empty condition at a next
time increment (e.g., output round) of the schedule clock. As
described above, a new finish time is computed for each respective
queue, and thus for each flow, when a new packet is received by the
respective queue. In this manner, the series of buckets are
progressively "emptied" as the schedule clock progresses, and new
buckets are filled as new queues receive new packets for
transmission and new associated empty times, and the buckets
progress from right to left as depicted in FIG. 6A.
Referring still to FIG. 6A, when the schedule clock advances to the
finish time of a bucket, the flows within the bucket, and thus
their queues, are regarded as completely serviced, and therefore,
empty. Those flows are considered inactive with respect to the link
when their queues are empty.
The schedule clock advances by one every time interval, T_Sclk,
given below:
##EQU00009## The schedule clock (represented as SCLK) advances
independently based on the flow activity on the corresponding link.
It should be noted that the SCLK does not necessarily advance at a
constant rate as a conventional clock.
.SIGMA.r.sub.i+CRR*.SIGMA.w.sub.I divided by the link capacity C,
represents the percentage of link usage at current CRR values. The
higher value of
.times. ##EQU00010## the slower the active SCLK advances. The
difference between the finish time of a queue and the schedule
clock represents the degree of backlog of the queue in terms of the
amount of time to empty the queue (empty time).
T.sub.empty=T.sub.finish-T.sub.SCLK
In addition to determining whether flows are active or inactive,
the schedule clock can also be used to pace flows to determine
whether any of them have exceeded their respective allocated
bandwidths. This can be done by ensuring T.sub.empty does not get
too large.
Referring still to FIG. 6A, being able to quickly determine which
VQs become empty simplifies the calculation of
.times..times..times..times..times..times..times. ##EQU00011## This
is due to the fact that the calculation can be done incrementally.
Given the old values, the new values can be calculated in one step
as follows:
.times..times..times..times..times..times. ##EQU00012##
.times..times..times..times. ##EQU00012.2##
.times. ##EQU00013## is also a term that can be computed
incrementally. When a flow moves from one bucket to another, its
r.sub.i and w.sub.i are subtracted from the sums of the old bucket
and added to that on the new buckets. For a flow that comes back
from a previously inactive (empty) state, its r.sub.i and w.sub.i
should be add to
.times..times..times..times..times..times..times. ##EQU00014##
too.
FIG. 6B graphically depicts the summation of all ri and wi in
accordance with one embodiment of the present invention. As shown
in FIG. 6B, the vertical axis is bandwidth and the horizontal axis
is time. The link capacity is as shown. The trace shows the
utilization of the link capacity as it changes over time (e.g., as
some flows become active and other flows become inactive).
Referring now to FIG. 7, a diagram of a bucket information base
(BIB) 700 in accordance with one embodiment of the present
invention is shown. The buckets depicted in FIG. 6A are implemented
as a series of counters within a database, the BIB 700, maintained
within each MPS. As depicted in FIG. 7, each bucket is implemented
as a ring total bandwidth counter and a corresponding ring total
weight counter. The counters are incremented to reflect the number
of flows, and their associated weights, within the bucket. The
schedule clock functions as a pointer that cycle through the
counters in accordance with the time increment at which their
respective flows will be empty, in the manner described above.
Thus, for example, at the next time increment, the schedule time
pointer will move to indicate the counters associated with bucket
1, and so on. In a preferred embodiment, BIB 700 is organized as a
two column and 8K long table as shown FIG. 7. BIB 700 is able to
sustain 16 accesses for every 50 ns, thereby allowing updates when,
for example, new packets arrive within the queues.
FIG. 8 shows a flow information base (FIB) 800 in accordance with
one embodiment of the present invention. An MPS uses the FIB 800 as
a flow descriptor. The FIB 800 contains fields specifying various
actions to be applied to the packets belonging to each flow
(transit forward on the ring, exit the ring to a specific port,
etc.) and fields holding flow parameters, such as r.sub.i and
w.sub.i. The finish time of a flow, which tracks its virtual queue
depth, is stored in the FIB. When packets arrive, the finish time
in the FIB is updated, and used to access the BIB as described
above. Thus the FIB is only accessed as packets arrive.
FIG. 9 shows a flow chart of the steps of an operating process 900
in accordance with one embodiment of the present invention. As
depicted in FIG. 9, process 900 shows the operating steps of an MPS
maintaining an accurate total of the amount of allocated bandwidth
on the network, as implemented within an MPTR.
Process 900 begins in step 901, where data packets for transmission
are received from a plurality of users by the queues of an MPS.
Within the MPS, the plurality of incoming packets from the various
users are assigned to a respective plurality of queues of the
MPS.
In step 902, data from the queues is routed onto the ring. Using a
fair arbitration scheme (e.g., round robin, etc.), a controller is
configured to empty the respective queues at a specified output
rate.
In step 903, a finish time is computed for each respect queue. The
finish time describes the time at which the respective queue will
be emptied using the current output rate.
In step 904, the queues are grouped into respective buckets based
on their respective finish times. To facilitate high-speed
tracking, the plurality of queues are grouped into multiple
buckets, or groups, in accordance with their respective finish
times. These groups are referred to as "buckets" due to the fact
that they include those queues having the same finish times. As
described above, the buckets can be implemented using respective
counter pairs within a database, the counter pairs configured to
track the total reserved r.sub.i having the same finish times and
their respective weights.
In step 905, a schedule clock is incremented in accordance with the
cycle time of the controller. As described above, a higher number
of active flows leads to a slower increment rate of the schedule
clock, and vice versa. The finish times are indexed with respect to
the schedule clock. The earliest bucket includes those queues
having a finish time indicating an empty condition at a first time
increment, the second earliest bucket includes those queues having
a finish time indicating an empty condition at a second time
increment later than the first time increment, and so on.
In step 906, the total r.sub.i of flows becoming inactive and their
associated weight are determined using the buckets. As described
above, counter pairs configured to track the reserved bandwidth of
queues having the same finish times and their respective weights
can be used to determine the allocated bandwidth of flows and their
associated weights becoming inactive on the next schedule clock
increment.
In step 907, determine the amount of unallocated bandwidth based
upon information obtained in step 906. As described above, the
amount of allocated bandwidth on the network is determined by
counting
.times..times..times..times..times..times..times. ##EQU00015## This
information allows the MPS to accurately determine an amount of
unallocated bandwidth available for distribution to the active
flows.
In step 908, new finish times are computed for the active flows as
new data arrives at the queues for transmission. Subsequently, in
step 909, process 900 continues by repeating steps 904-909. In this
manner, the series of buckets are progressively "emptied" as the
schedule clock progresses, and new buckets are filled as new queues
receive new packets for transmission and new associated empty
times.
Thus, the determination of the amount of allocated bandwidth can be
accomplished in real time, thereby allowing the efficient
allocation of unallocated bandwidth in real time while maintaining
quality of service. The earliest bucket (e.g., bucket 0) shows all
queues which will be empty in the next time increment. In so doing,
the present invention enables the efficient allocation of available
bandwidth, since the MPS is capable of tracking total allocated
bandwidth in real time. This allows the efficient allocation of
unused bandwidth in real time while maintaining QoS.
The foregoing descriptions of specific embodiments of the present
invention have been presented for purposes of illustration and
description. They are not intended to be exhaustive or to limit the
invention to the precise forms disclosed, and obviously many
modifications and variations are possible in light of the above
teaching. The embodiments were chosen and described in order to
best explain the principles of the invention and its practical
application. It is intended that the scope of the invention be
defined by the Claims appended hereto and their equivalents.
* * * * *